Methods and systems for real-time continuous updates

ABSTRACT

Embodiments of the present invention provide fine grain concurrency control for transactions in the presence of database updates. During operations, each transaction is assigned a snapshot version number or SVN. A SVN refers to a historical snapshot of the database that can be created periodically or on demand. Transactions are thus tied to a particular SVN, such as, when the transaction was created. Queries belonging to the transactions can access data that is consistent as of a point in time, for example, corresponding to the latest SVN when the transaction was created. At various times, data from the database stored in a memory can be updated using the snapshot data corresponding to a SVN. When a transaction is committed, a snapshot of the database with a new SVN is created based on the data modified by the transaction and the snapshot is synchronized to the memory. When a transaction query requires data from a version of the database corresponding to a SVN, the data in the memory may be synchronized with the snapshot data corresponding to that SVN.

CROSS-REFERENCE TO RELATED APPLICATIONS

This application is a divisional application of U.S. patent application Ser. No. 12/144,486; entitled “METHODS AND SYSTEMS FOR REAL-TIME CONTINUOUS UPDATES” by Kapil Surlaker, Ravi Krishnamurthy, Krishnan Meiyyappan, Alan Beck, Hung Tran, Jeremy Branscome, and Joseph Chamdani; filed on Jun. 23, 2008.

This application is related to the following U.S. Patent Applications and Patents, which are herein incorporated by reference in their entirety: U.S. patent application Ser. No. 11/895,952, filed on Aug. 27, 2007, entitled Methods and Systems for Hardware Acceleration of Database Operations and Queries, by Joseph I. Chamdani et al.; U.S. patent application Ser. No. 11/895,998, filed on Aug. 27, 2007, entitled Hardware Acceleration Reconfigurable Processor for Accelerating Database Operations and Queries, by Jeremy Branscome et al.; U.S. patent application Ser. No. 11/895,997, filed on Aug. 27, 2007, entitled Processing Elements of a Hardware Acceleration Reconfigurable Processor for Accelerating Database Operations and Queries, by Jeremy Branscome et al.; U.S. patent application Ser. No. 12/168,821, filed on Jul. 7, 2008, entitled Methods and Systems for Generating Query Plans that are Compatible for Execution in Hardware, by Ravi Krishnamurthy et al.; U.S. patent application Ser. No. 12/099,076, filed on Apr. 7, 2008, entitled Methods and Systems for Run-Time Scheduling Database Operations that are Executed in Hardware, by Joseph I. Chamdani et al.; U.S. patent application Ser. No. 12/099,131, filed on Apr. 7, 2008, entitled Accessing Data in a Column Store Database Based on Hardware Compatible Data Structures, by Liuxi Yang et al.; U.S. patent application Ser. No. 12/099,133, filed on Apr. 7, 2008, entitled Accessing Data in a Column Store Database Based on Hardware Compatible Indexing and Replicated Reordered Columns, by Krishnan Meiyyappan et al.; and U.S. patent application Ser. No. 12/144,303, filed on Jun. 23, 2008, entitled Fast Bulk Loading and Incremental Loading of Data into a Database, by James Shau et al.

BACKGROUND

Despite their different uses, applications, and workload characteristics, most systems run on a common Database Management System (DBMS) using a standard database programming language, such as Structured Query Language (SQL). Most modern DBMS implementations (Oracle, IBM DB2, Microsoft SQL, Sybase, MySQL, PostgreSQL, Ingress, etc.) are implemented on relational databases, which are well known to those skilled in the art.

Typically, a DBMS has a client side where applications or users submit their queries and a server side that executes the queries. On the server side, most enterprises employ one or more general purpose servers. However, although these platforms are flexible, general purpose servers are not optimized for many enterprise database applications. In a general purpose database server, all SQL queries and transactions are eventually mapped to low level software instructions called assembly instructions, which are then executed on a general purpose microprocessor (CPU). The CPU executes the instructions, and its logic is busy as long as the operand data are available, either in the register file or on-chip cache. To extract more parallelism from the assembly code and keep the CPU pipeline busy, known CPUs attempt to predict ahead the outcome of branch instructions and execute down the SQL code path speculatively. Execution time is reduced if the speculation is correct; the success of this speculation, however, is data dependent. Other state-of-the-art CPUs attempt to increase performance by employing simultaneous multithreading (SMT) and/or multi-core chip multiprocessing (CMP). To take advantage of these, changes have to be made at the application or DBMS source code to manually create the process/thread parallelism for the SMT or CMP CPUs. This is generally considered highly undesirable.

Unfortunately, general purpose CPUs are not efficient for database applications. Branch prediction is generally not accurate because database processing involves tree traversing and link list or pointer chasing that is very data dependent. Known CPUs employ the well-known code-flow (or Von Neumann) architecture, which uses a highly pipelined instruction flow (rather than a dataflow where operand data is pipelined) to operate on data stored in the CPUs tiny register files. Real database workloads, however, typically require processing Gigabytes to Terabytes of data, which overwhelms these tiny registers with loads and reloads. On-chip cache of a general purpose CPU is not effective since it's relatively too small for real database workloads. This requires that the database server frequently retrieve data from its small memory or disk. Accordingly, known database servers rely heavily on squeezing the utilization of their small system memory size and disk input/output (I/O) bandwidth. Those skilled in the art recognize that these bottlenecks between storage I/O, the CPU, and memory are very significant performance factors.

However, overcoming these bottlenecks is a complex task because typical database systems consist of several layers of hardware, software, etc., that influence the overall performance of the system. These layers comprise, for example, the application software, the DBMS software, operating system (OS), server processor systems, such as its CPU, memory, and disk I/O and infrastructure. Traditionally, performance has been optimized in a database system horizontally, i.e., within a particular layer. For example, many solutions attempt to optimize various solutions for the DBMS query processing, caching, the disk I/O, etc. These solutions employ a generic, narrow approach that still fails to truly optimize the large performance potentials of the database system, especially for relational database systems having complex read-intensive applications. In addition, in order to support updates, the conventional solutions typically employ complex mechanisms to provide versioning control and concurrent access.

BRIEF DESCRIPTION OF THE DRAWINGS

The accompanying drawings, which are incorporated in and constitute a part of this specification, illustrate embodiments of the invention and together with the description, serve to explain the principles of the invention. In the figures:

FIG. 1 illustrates an exemplary system that is consistent with the principles of the present invention;

FIG. 2 illustrates exemplary system topologies that are consistent with the principles of the present invention;

FIG. 3A illustrates a database system and FIG. 3B illustrates some of the optimizations of the present disclosure;

FIG. 4 illustrates a functional architecture of the custom computing (C2) software of the present invention;

FIG. 5 illustrates a protocol stack employed by the C2 software and a Hardware Acceleration Reconfigurable Processor (HARP) of the present invention;

FIG. 6 illustrates an exemplary architecture of a HARP;

FIG. 7 illustrates a column store database and associated data structures employed by some embodiments of the present invention;

FIG. 8 illustrates a table column layout and associated data structures employed by some embodiments of the present invention;

FIG. 9 illustrates an exemplary machine code database instruction flow for a SQL query that is consistent with the principles of the present invention;

FIG. 10 illustrates an exemplary dataflow for a SQL query through processing elements in the HARP in accordance with the principles of the present invention;

FIG. 11 illustrates an exemplary logic flow of the addressing schemes employed in the present invention;

FIG. 12 illustrates a structure of the global database virtual address scheme of the present invention;

FIG. 13 illustrates an exemplary timeline of transactions and snapshots to support multi-versioning and continuous updates;

FIG. 14 illustrates exemplary synching of pages between different snapshots;

FIG. 15 illustrates an example of synching a page from one snapshot to a subsequent snapshot;

FIG. 16 illustrates an example of synching a page from one snapshot to a previous snapshot; and

FIGS. 17 and 18 illustrate associative accumulation machine code database instructions that may be employed to change page contents from one snapshot version to another.

DETAILED DESCRIPTION

Embodiments of the present invention provide fine grain concurrency control for transactions in the presence of database updates. During operations, each transaction is assigned a snapshot version number or SVN. A SVN refers to a historical snapshot of the database that can be created periodically or on demand. Transactions are thus tied to a particular SVN, such as, when the transaction was created. Queries belonging to the transactions can access data that is consistent as of a point in time, for example, corresponding to the latest SVN when the transaction was created.

At various times, data from the database stored in a memory can be updated using the snapshot data corresponding to a SVN. For example, when a transaction is committed, a snapshot of the database with a new SVN is created based on the data modified by the transaction and the snapshot is synchronized to the memory. As another example, when a transaction query requires data from a version of the database corresponding to a SVN, the data in the memory may be synchronized with the snapshot data corresponding to that SVN. This feature, in essence, creates multiple versions of the database and allows read-intensive database applications to operate in the presence of database updates.

Due to the comprehensive nature of the present inventions in the C2 solution, the figures are presented generally from a high level of detail and progress to a low level of detail. For example, FIGS. 1-3 illustrate exemplary systems and topologies enabled by the present invention. FIG. 4-5 illustrate the architecture of the C2 software. FIG. 6 illustrates the architecture of a HARP module. FIGS. 7-8 illustrate the database format and data structures employed by the C2 solution of the present invention. FIGS. 9-10 illustrates an example execution of a SQL query by the C2 solution of the present invention. Reference will now be made in detail to the exemplary embodiments of the invention, which are illustrated in the accompanying drawings. FIG. 11 illustrates an exemplary logic flow of the addressing schemes employed. FIG. 12 illustrates a structure of the global database virtual address scheme. FIG. 13 illustrates an exemplary timeline of transactions and snapshots to support multi-versioning and continuous updates. FIG. 14 illustrates exemplary synching of pages between different snapshots. FIG. 15 illustrates an example of synching a page from one snapshot to a subsequent snapshot. FIG. 16 illustrates an example of synching a page from one snapshot to a previous snapshot. FIGS. 17 and 18 illustrate associative accumulation machine code database instructions that may be employed to change page contents from one snapshot version to another. Wherever possible, the same reference numbers will be used throughout the drawings to refer to the same or like parts.

FIG. 1—An Exemplary C2 System

The present invention employs a custom computing (C2) solution that provides a significant gain in performance for enterprise database applications. In the C2 solution, a node or appliance may comprise the host (or base) system that is combined with hardware acceleration reconfigurable processors (HARP). These HARPs are specially designed to optimize the performance of database systems and its applications, especially relational database systems and read-intensive applications.

A host system may be any standard or pre-existing DBMS system. In general, such systems will comprise a standard general purpose CPU, a system memory, I/O interfaces, etc.

The HARPs are coupled to the host system and are designed to offload repetitive database operations from the DBMS running on the host system. The HARPs utilize dataflow architecture processing elements that execute machine code instructions that are defined for various database operations. The C2 solution may employ a node that is scalable to include one HARP, or multiple HARPs. In addition, the C2 solution may use a federated architecture comprising multiple nodes, i.e., multiple DBMS servers that are enhanced with the C2 solution.

In some embodiments, the C2 solution employs an open architecture and co-processor approach so that the C2 hardware can be easily integrated into existing database systems. Of note, the hardware acceleration of the C2 solution utilizes novel machine code database instructions to execute certain fragments of a query in a dataflow and using parallel, pipelined execution.

In the present invention, the C2 solution also comprises software that orchestrates the operations of the DBMS running on the host system and the HARPs. The C2 software is configured with a flexible, layered architecture to make it hardware and database system agnostic. Thus, the C2 software is capable of seamlessly working with existing DBMSs based on this open architecture.

In general, the C2 software receives the query from the DBMS and breaks the query down into query fragments. The C2 software then decides which of these query fragments can be appropriately handled in software (in the C2 software itself or back in the originating DBMS) or, ideally, with hardware acceleration in the HARPs. All or part of the query may be processed by the C2 software and HARPs.

In addition, in order to maximize the efficiency of the hardware acceleration, the C2 solution stores its databases in compressed, column-store format and utilizes various hardware-friendly data structures. The C2 solution may employ various compression techniques to minimize or reduce the storage footprint of its databases. The column-store format and hardware-friendly data structures allow the HARPs or C2 software to operate directly on the compressed data in the column-store database. The column-store database may employ columns and column groups that are arranged based on an implicit row identifier (RID) scheme and RID to primary column to allow for easy processing by the HARPs. The hardware-friendly data structures also allow for efficient indexing, data manipulation, etc. by the HARPs.

For example, the C2 solution utilizes a global virtual address space for the entire database to greatly simplify and maximize efficiency of create, read, update, and delete operations of data in a database. In some embodiments, the columns and column groups are configured with a fixed width to allow for arithmetic memory addressing and translation from a virtual address to a physical memory address. On-demand and speculative prefetching may also be utilized by the C2 solution to hide I/O bandwidth latency and maximize HARP utilization.

Referring now to FIG. 1, an exemplary system 1 f the C2 solution is illustrated. As shown, system 100 may comprise an application 102 that is running on a client 104, such as a personal computer or other system. Application 102 interfaces a DBMS 106 across a network 108, such as the Internet, local area network, etc. DBMS 106 may further interface one or more databases stored in storage infrastructure 112. For purposes of explanation, DBMS 106 and its components may be collectively referred to in this disclosure as a node of system 100. Although FIG. 1 shows a single node, system 100 may of course comprise multiple nodes. The various components of FIG. 1 will now be further described.

Application 102 may be any computer software that requests the services of DBMS 106. Such applications are well known to those skilled in the art. For example, application 102 may be a web browser in which a user is submitting various search requests. Of course, application 102 may be another system or software that is consuming the services of DBMS 106 and submitting queries to DBMS 106.

Client 104 represents the hardware and software that supports the execution of application 102. Such clients are well known to those skilled in the art. For example, client 104 may be a personal computer or another server.

DBMS 106 is any computer software that manages databases. In general, DBMS 106 controls the organization, storage, management, retrieval of data in a database. As is well known, these types of systems are common for supporting various SQL queries on relational databases (and thus may also be known as a RDBMS). Due to its open architecture, various DBMS systems may be employed by the present invention. Typical examples of DBMSs include Oracle, DB2, Microsoft Access, Microsoft SQL Server, PostgreSQL, and MySQL.

In some embodiments, and for purposes of explanation, DBMS 106 is shown comprising C2 software 110 interfacing MySQL software 114 via an API 116. MySQL software 114 is open source software that is sponsored and provided by MySQL AB and is well known to those skilled in the art. Of course, any DBMS software, such as those noted above, may be employed in the present invention.

C2 software 110 orchestrates the execution of a query forwarded from DBMS 106, and thus, operates in conjunction with MySQL software 114. For example, in the C2 software 110, SQL queries are broken down into query fragments and then routed to the most appropriate resource. A query fragment may be handled in C2 hardware, i.e., HARP module 204. (HARP module 204 is further described with reference to FIG. 2.) The query fragment may also be processed in the C2 software itself, or returned for handling by MySQL software 114.

In general, C2 software 110 utilizes a flexible, layered architecture to make it hardware and database system agnostic. For example, C2 software 110 may operate as a storage engine of MySQL software 114. As is well known, MySQL software 114 may provide an API 116 for storage engines, which C2 software 110 may plug in to. API 116 comprises the software that specifies how the C2 software 110 and MySQL software 114 will interact, how they will request services from each other, such as SQL queries and results.

As a storage engine, C2 software 110 may employ the MySQL API 116 to provide various storage mechanisms, indexing facilities, locking levels, and ultimately provide a range of different functions and capabilities that are transparent to MySQL software 114. As noted above, this is one aspect of how the present invention overcomes the generic approach in known solutions without having to sacrifice performance for functionality, or fine tune the database. Of note, although FIG. 1 shows a single storage engine, MySQL software 114 may be coupled to multiple storage engines (not shown) in addition to C2 software 110. C2 software 110 is also described in further detail with reference to FIGS. 4-5.

Network 108 represents the communication infrastructure that couples application 102 and DBMS 106. For example, network 108 may be the Internet. Of course, any network, such as a local area network, wide area network, etc., may be employed by the present invention.

Storage infrastructure 112 comprises the computer storage devices, such as disk arrays, tape libraries, and optical drives that serve as the storage for the databases of system 100. Storage infrastructure 112 may employ various architectures, such as a storage area network, network attached storage, etc., which are known to those skilled in the art.

In some embodiments, the C2 solution stores its databases in storage infrastructure 112 in column-store format. Column-store format is where data is stored in columns or groups of columns. Column-store format is advantageous for data fetching, scanning, searching, and data compression. The column-store format may employ fixed width columns and column groups with implicit RIDs and a RID to primary key column to allow for arithmetic memory addressing and translation. This allows HARPs 204 to utilize hardware processing for database processing, such as column hopping, and to operate directly on the compressed data in the columns.

In contrast, in typical DBMS environments, data is stored in row-store format. Row-store format is sometimes considered by those skilled in the art for having better performance in data updates and record retrieval; thus, it is sometimes considered to have better functionality over column-store databases in most applications with a high ratio of updates over reads. In the present invention, however, the C2 solution achieves better performance by using hardware acceleration with a column-store database, yet it still delivers the functionality and benefits of row-store databases. The column store format used by the C2 solution of the present invention is further described with reference to FIGS. 7-8.

FIG. 2—System Topologies

FIG. 2 illustrates exemplary system topologies that are consistent with the principles of the present invention. As shown, FIG. 2 illustrates a basic C2 node topology, a scale up C2 node topology, and a scale out topology. These various topologies may be utilized to customize the C2 solution for various sizes of databases and desired performance. In addition, these topologies are provided to illustrate that the C2 solution can be easily scaled up to virtually any size of database or performance.

First, the basic C2 node will be explained, which comprises a single host system 202 and a single HARP module 204. Variations of this basic node will then be explained to show how the basic node can be scaled up and how multiple nodes can be employed in a federated architecture.

The basic C2 node topology may comprise a host system 202 and a hardware acceleration reconfigurable processor (HARP) module 204. Collectively, host 202 and HARP module 204 may be referred to as a node or appliance. In some embodiments, host system 202 and HARP module 204 are coupled together over a known communications interface, such as a PCIe or hypertransport (HT) interface. In terms of packaging, host system 202 and HARP module 204 may be built on one or more cards or blades that are bundled together in a common chassis or merely wired together. In the C2 solution, host system 202 and HARP module 204 may be flexibly packaged using a modular form factor for ease of installation and scaling.

The host system 202 may comprise a general purpose CPU, such as a Xeon x86 processor by the Intel Corporation, and a memory, such as a dynamic random access memory. Such types of host systems are well known to those skilled in the art. In general, in the C2 solution, host system 202 will be used to process parts of a query that are less time consuming (i.e., slow path portion), such as server-client connection, authentication, SQL parsing, logging, etc. However, in order to optimize performance, the bulk of query execution (i.e., the fast path portion) is offloaded to the HARP module 204.

Host system 202 may run MySQL software 114 and also run C2 software 110 that orchestrates query processing between MySQL 114 and HARP 204. In particular, C2 software 110 will decompose a query into a set of query fragments. Each fragment comprises various tasks, which may have certain dependencies. C2 software 110 will determine which fragments and tasks are part of the fast path portion and offload them to the HARP module 204. Appropriate tasks for the selected query fragments are sent to HARP module 204 with information on the database operation dependency graph. Within the HARP module 204, tasks are further broken down into parallel/pipelined machine code operations (known as MOPs) and executed in hardware.

HARP module 204 comprises processing logic (HARP logic 302) and a relatively large memory (HARP memory 304) for hardware accelerating database operations of the node. In some embodiments, HARP module 204 is configured to handle various repetitive database tasks, such as table scanning, indexing, etc. In the C2 solution, HARP module 204 can receive high-level database query tasks (not just low-level read/write or primitive computation tasks as is typically for a general purpose processor) in the form of machine code database instructions.

HARP logic 302 is the hardware that executes machine code database instructions for the database tasks being handled by HARP module 204. To adapt to application requirement changes, the HARP logic 302 is designed to have hardware re-configurability. Accordingly, in some embodiments, HARP logic 302 is implemented using field programmable gate arrays (FPGAs). However, any type of custom integrated circuit, such as application specific integrated circuits (ASICs), may be implemented as HARP logic 302.

HARP memory 304 serves as the memory of HARP module 204. In order to maximize the efficiency of the HARP logic 302, the HARP memory 304 may be implemented using relatively large amounts of memory. For example, in some embodiments, the HARP memory 304 in a HARP module 204 may comprise 256 Giga-Bytes or more of RAM or DRAM. Of course, even larger amounts of memory may be installed in HARP module 204. HARP logic 302 and HARP memory 304 are further described with reference to FIG. 6.

In addition to the basic C2 node, a scale up C2 node topology may be used as an extension of the basic C2 node. As shown, host system 202 may now be coupled to a plurality or array of 1-N HARP modules 204. In this type of node, a PCIe switch or other suitable switching fabric may couple these components together with storage infrastructure 112. Of course, other internal arrangements for a scale up C2 node may be utilized in the present invention.

Going further, a scale out topology can be used for multiple C2 nodes. As shown, the scale out topology may comprise various combinations of either the basic or scale up C2 nodes. For example, as shown, the scale out topology may comprise Nodes 1-M, which are coupled to storage infrastructure 112. In FIG. 2, Node 1 is shown as a basic C2 node, while Node M is shown as a scale up node. A control node 206 is also shown and manages the operations of Nodes 1-M. Control node 206 is shown as a separate node; however, those skilled in the art will recognize the role of control node 206 by any of Nodes 1-M. Other variations in node hierarchy and management are within the scope of the present invention. Of course, this topology may also comprise a variety of combinations of nodes.

FIGS. 3A and 3B—Some Advantages of the Present Invention

FIG. 3A illustrates a prior art database system and FIG. 3B illustrates an exemplary implementation of the C2 solution for the present invention. In FIG. 3A, a typical prior art database system is shown. An SQL query is submitted to a DBMS (e.g., MySQL), which runs on top of a typical operating system. The CPU attempts to then execute the SQL query. However, because the CPU is a general purpose CPU it executes this query based on software, which has several limitations.

In contrast, as shown in FIG. 3B, the SQL query may submitted to a C2 system having a DBMS that comprises a top layer DBMS software (i.e., MySQL) 114 and C2 software 110. C2 software 110 interfaces with the DBMS software 114 to orchestrate and optimize processing of the SQL query.

In particular, C2 software 110 may identify portions of the query, i.e., the fast path portion, which is better handled in hardware, such as HARP module 204. Such portions may be those fragments of the query that are repetitive in nature, such as scanning, indexing, etc. In the prior art system, the DBMS is limited by its own programming, the operating system, and the general purpose CPU. The present invention avoids these bottlenecks by offloading fast path portions of a query to HARP module 204.

As shown, HARP module 204 comprises HARP logic 302 and a HARP memory 304 to accelerate the processing of SQL queries. In order maximize the use of HARP module 204, the present invention may also utilize column store databases. Whereas the prior art system is hindered by the limitations of a standard row store database. These features also allow the present invention to maximize the performance of the I/O between the operating system and storage.

For ease of implementation, C2 software 110 may be implemented on well-known operating systems. The operating system will continue to be used to perform basic tasks such as controlling and allocating memory, prioritizing system requests, controlling input and output devices, facilitating networking, and managing files and data in storage infrastructure 112. In some embodiments, various operating systems, such as Linux, UNIX, and Microsoft Windows, may be implemented.

FIGS. 3A and 3B are provided to illustrate some of the differences between the present invention and the prior art and advantages of the present invention. Those skilled in the art will also recognize that other advantages and benefits may be achieved by the embodiments of the present invention. For purposes of explanation, the present disclosure will now describe the C2 software, hardware, data structures, and some operations in further detail.

FIG. 4—C2 Software Architecture

As noted, C2 software 110 orchestrates the processing a query between MySQL software 114 and HARP module 204. In some embodiments, C2 software 110 runs as an application on host system 202 and as a storage engine of MySQL software 114. FIG. 4 illustrates an architecture of the C2 software 110. As shown, C2 software 110 comprises a query and plan manager 402, a query reduction/rewrite module 404, an optimizer 406, a post optimizer module 408, a query plan generator 410, an execution engine 412, a buffer manager 414, a task manager 416, a memory manager 418, a storage manager 420, an answer manager 422, an update manager 424, shared utilities 426, and a HARP manager 428. Each of these components will now be briefly described.

Query and plan manager 402 analyzes and represents the query received from the MySQL software 114, annotates the query, and provides a representation of the query plan. Query reduction/rewrite module 404 breaks the query into query fragments and rewrites the query fragments into tasks. Rewrites may be needed for compressed domain rewrites and machine code database instruction operator rewrites. Optimizer 406 performs cost-based optimization to be done using cost model of resources available to C2 software 110, i.e., HARP module 204, resources of C2 software 110 itself using software operations, or MySQL software 114.

These modules interact with each other to determine how to execute a query, such as a SQL query from MySQL software 114. The data structures output by the query plan generator 410 will be the same data structure that the optimizer 406 and the rewrite module 404 will operate on. Once a parsed SQL query has been represented in this data structure (converted, for example, from MySQL), manager 402 rewrites the query such that each fragment of the query can be done entirely in MySQL software 114, in C2 software 110, or in HARP module 204. Once the final query representation is available, the rewrite module 404 goes through and breaks the graph into query fragments.

Post optimizer module 408 is an optional component that rewrites after the optimizer 406 for coalescing improvements found by optimizer 406. Query plan generator 410 generates an annotations-based, template-driven plan generation for the query tasks. Execution engine 412 executes the query fragments that are to be handled by software or supervises the query execution in HARP module 204 via HARP manager 428.

Buffer manager 414 manages the buffers of data held in the memory of host 202 and for the software execution tasks handled by host 202. Task manager 416 orchestrates the execution of all the tasks in HARP module 204 and software, i.e., in execution engine 412 or MySQL software 114.

Memory manager 416 manages the virtual address and physical address space employed by C2 software 110 and HARP module 204 in HARP memory 304. In some embodiments, memory manager 416 utilizes a 50 bit VA addressing (i.e., in excess of 1 petabyte). This allows C2 software 110 to globally address an entire database and optimize hardware execution of the query tasks.

Storage manager 420 is responsible for managing transfers of data from HARP memory 304 to/from storage infrastructure 112. Answer manager 422 is responsible for compiling the results of the query fragments and providing the result to MySQL software 114 via the API 116.

Update manager 424 is responsible for updating any data in the database stored in storage infrastructure 112. Shared utilities 426 provide various utilities for the components of C2 software 110. For example, these shared utilities may include a performance monitor, a metadata manager, an exception handler, a compression library, a logging and recovery manager, and a data loader.

HARP manager 428 controls execution of the tasks in HARP module 204 by setting up the machine code database instructions and handles all interrupts from any of the hardware in HARP module 204. In some embodiments, HARP manager 428 employs a function library known as a Hardware Acceleration Function Library (HAFL) in order to make its function calls to HARP module 204.

FIG. 5—Protocol Stack of C2 Software

As shown, a SQL query is received in the RDBMS layer, i.e., MySQL software 114. MySQL software 114 then passes the SQL query via API 116 to C2 software 110. In C2 software 110, the SQL query is processed and executed. At this layer, C2 software 110 also manages retrieving data for the SQL query, if necessary, from storage infrastructure 112 or from host system 202.

In order to communicate with HARP module 204, HARP manager 428 employs the HAFL layer in order to make its function calls to HARP module 204. In order to allow for variances in hardware that may exist in HARP module 204, the protocol stack may also comprise a hardware abstraction layer. Information is then passed from C2 software 110 to HARP module 204 in the form of machine code database instructions via an interconnect layer. As noted, this interconnect layer may be in accordance with the well-known PCIe or HT standards.

Within HARP module 204, the machine code database instructions are parsed and forwarded to HARP logic 302. These instructions may relate to a variety of tasks and operations. For example, as shown, the protocol stack provides for systems management, task coordination, and direct memory access to HARP memory 304. In HARP logic 302, machine code database instructions are interpreted for the various types of processing elements (PE). HARP logic 302 may interface with HARP memory 304, i.e., direct memory access by utilizing the memory management layer.

FIG. 6—HARP Logic

FIG. 6 illustrates an exemplary architecture of the HARP logic 302. As shown, HARP logic 302 may comprise a set of processing cores 602, 604, 606, and 608, and switching fabric 610. Processing core 602 (as well as cores 604, 606, and 608) may comprise a set of processing elements (PEs) 620. In the embodiment shown, processing cores 602, 604, 606, and 608 each comprise two PEs; of course, each processing core may comprise any number of PEs.

In addition to its PEs, processing core 602 may comprise a task processor 612, a memory manager 614, a buffer cache 616, and an interconnect 618. One or more these components may be duplicated or removed from the other processing cores 604, 606, and 608. For example, as shown, core 602 may be the sole core that includes task processor 612 and an interconnect 618. This architecture may be employed because cores 602, 604, 606, and 608 are connected via switching fabric 610 and may operate logically as a single processor or processor core. Of course, one skilled in the art will recognize that various redundancies may be employed in these processing cores as desired.

Task processor 612 is the hardware that supervises the operations of the processing cores 602, 604, 606, and 608. Task Processor 612 is a master scheduling and control processing element, disconnected from the direct dataflow of the execution process for a query. Task processor 612 maintains a running schedule of machine code database instructions which have completed, are in progress, or are yet to execute, and their accompanying dependencies, the Task Processor 612 may also dispatch machine code database instructions for execution and monitor their progress. Dependencies can be implicit, or explicit in terms of strong intra- or inter-processor release criteria. Machine code database instructions stalled for software-assist can be context-switched by the Task Processor 612, which can begin or continue execution of other independent query tasks, to optimize utilization of execution resources in HARP logic 302.

Memory manager 614 is the hardware that interfaces HARP memory 304. For example, memory manager 614 may employ well known memory addressing techniques, such as translation look-aside buffers to map the global database virtual address space to a physical address in HARP memory 304 to access data stored in HARP memory 304.

Buffer cache 616 serves as a small cache for a processing core. For example, temporary results or other meta-data may be held in buffer cache 616.

PCIe interconnect 618 is the hardware that interfaces with host system 202. As noted, interconnect 618 may be a PCIe or HT interconnect.

PEs 620 represent units of the hardware and circuitry of HARP logic 302. As noted, PEs 620 utilize a novel dataflow architecture to accomplish the query processing requested of HARP logic 302. In particular, PEs 620 implement execution of an assortment of machine code database instructions that are known as Macro Ops (MOPs) and Micro Ops (UOPs). MOPs and UOPs are programmed and executed by the PEs 620 to realize some distinct phase of data processing needed to complete a query. MOPs and UOPs are just example embodiments of machine code database instructions; other types of instruction sets for high level database operations of course may be used by the C2 solution.

PEs 620 pass logical intermediate MOP results among one another through a variable-length dataflow of dataflow tokens, carried across an interconnect data structure (which is a physical data structure and not a software data structure) termed an Inter-Macro Op Communication (IMC) path. Of note, the IMC paths and self routing fabric 610 allow HARP module 204 to utilize a minimal amount of reads/writes to HARP memory 304 by keeping most intermediate results flowing through the IMCs in a pipelined, parallel fashion. IMC may be temporarily stored in buffer caches 616 and interconnect fabric 610; however, IMCs can also be dispatched out through interconnect 618 to other PEs 620 on another HARP module.

In the dataflow concept, each execution step, as implemented by a MOP and its accompanying UOP program, can apply symmetrically and independently to a prescribed tuple of input data to produce some tuple of result. Given the independence and symmetry, any number of these tuples may then be combined into a list, matrix, or more sophisticated structure to be propagated and executed in pipelined fashion, for optimal execution system throughput. These lists of tuples, comprised fundamentally of dataflow tokens, are the intermediate and final results passed dynamically among the MOPs via IMC.

Although the dataflow travels over physical links of potentially fixed dimension, the logical structure of the contents can be multi-dimensional, produced and interpreted in one of two different ways: either with or without inherent, internal formatting information. Carrying explicit internal formatting information allows compression of otherwise extensive join relationships into nested sub list structures which can require less link bandwidth from fabric 610 and intermediate storage in buffer cache 616, at the cost of the extra formatting delimiters, increased interpretation complexity and the restriction of fixing the interpretation globally among all consumers. Without inherent formatting, a logical dataflow may be interpreted by the consumer as any n-dimensional structure having an arbitrary but consistent number of columns of arbitrary but consistent length and width. It should be noted that the non-formatted form can be beneficial not only in its structural simplicity, but in the freedom with which consumer MOPs may interpret, or reinterpret, its contents depending upon the purpose of the execution step a consumer is implementing.

The dataflow used in realizing a given query execution can be described by a directed acyclic graph (DAG) with one intervening MOP at each point of flow convergence and bifurcation, one MOP at each starting and ending point, as well as any point necessary in between (i.e. single input & output MOP). The DAG must have at least one starting and one ending point, although any larger number may be necessary to realize a query. MOPs which serve as the starting point are designed to begin the dataflow by consuming and processing large amounts of data from local storage. Ending point MOPs may terminate the dataflow back into local storage, or to a link which deposits the collected dataflow (result table list) into host CPU memory. An example of a DAG for a well-known TPC-H query is shown in FIG. 9.

As mentioned above, MOP DAGs can physically and logically converge or bifurcate, programmatically. The physical convergence is accomplished with a multi-input MOPs which relate inputs in some logical fashion to produce an output comprised of all inputs (e.g. composition, merge, etc.). The physical bifurcation is accomplished by means of multicast technology in the IMC fabric, which dynamically copies an intermediate result list to multiple consumer MOPs. These mechanisms work together to allow realization of any desired DAG of MOP execution flow.

In the present invention, each MOP is configured to operate directly on the compressed data in the column-store database and realizes some fundamental step in query processing. MOPs are physically implemented and executed by PEs 620 which, depending on specific type, will realize a distinct subset of all MOP types. MOPs work systematically on individual tuples extracted either from local database storage in HARP memory 304 or the IMC dataflow, producing output tuples which may be interpreted by one or more MOP processes downstream.

UOPs are the low-level data manipulators which may be combined into a MOP-specific UOP program accompanying a MOP, to perform analysis and/or transformation of each tuple the MOP extracts. MOPs which utilize UOP programs are aware of the dependency, distributing selected portions of each tuple to the underlying UOP engine, extant within all PEs 620 supporting such MOPs. For each set of inputs from each tuple, the UOP program produces a set of outputs, which the MOP may use in various ways to realize its function.

For example, one manner a MOP may use UOP output is to evaluate each tuple of a list of tuples for a set of predicating conditions, where the MOP decides either to retain or to drop each tuple based on the UOP result. Another manner is for the UOP to perform an arithmetic transformation of each input tuple, where the MOP either appends the UOP result to form a larger logical tuple, or replaces some portion of the input tuple to form the output tuple.

Given a finite number of execution resources in PEs 620, the full MOP dataflow DAG needed to execute a query may be partitioned into segments of connected MOPs called tasks. These tasks are then scheduled by task processor 612 for execution in a sequential fashion, as MOP execution resources become available in PEs 620. Significant in this process is the propagation of the execution dataflow among these tasks, such that the entire query result is accurately and consistently computed, regardless of how each task is apportioned and regardless of the latency between scheduling each task.

One method that may be employed in HARP logic 302 is to treat each task atomically and independently, terminating the dataflow back into local storage in HARP memory 304 at the end of each task and restarting that dataflow at the beginning of the subsequent task by reloading it from HARP memory 304. In some embodiments, a more efficient method may be employed to pipeline tasks at their finer, constituent MOP granularity, where at least one MOP of a new task may begin execution before all MOPs of the previous task have finished. This fine-grained method is referred to as Task Pipelining.

Keeping the dataflow alive over task boundaries is a key to realizing the extra efficiency of Task Pipelining. To accomplish this in the C2 solution, IMCs may include the ability to dynamically spill, or send their dataflow to an elastic buffer backed by HARP memory 304, pending the awakening of a consumer MOP which will continue the dataflow. On scheduling the consumer MOP, IMCs are able to fill dynamically, reading from the elastic buffer in HARP memory 304 as necessary to continue execution, pulling out any slack that may have built up in the dataflow while waiting for the scheduling opportunity. Task Pipelining with these mechanisms then may provide a more efficient use of execution resources, down to the MOP granularity, such that a query may be processed as quickly as possible.

High-latency, low-bandwidth, non-volatile storage in storage infrastructure 112 often holds the contents of a query, due to the sheer volume of data involved. Because execution rates can outstrip the bandwidth available to read from such storage, tasks requiring latent data can shorten execution time by starting and progressing their dataflow execution at the rate the data arrives, instead of waiting for an entire prefetch to complete before beginning execution. This shortcut is referred to as Prefetch Pipelining. The C2 solution may employ both on-demand prefetching and speculative prefetching. On-demand prefetching is where data is prefetched based on the progress of the dataflow. Speculative prefetching is where data is prefetched based on an algorithm or heuristic that estimates the data is likely to be requested as part of a dataflow.

In the present invention, realizing Prefetch Pipelining is accomplished by having one or more MOPs beginning a task's dataflow are capable of accepting data progressively as it is read from slow storage in storage infrastructure 112. IMCs are capable of filling progressively as data arrives, as are all MOPs already designed to read from local storage in HARP memory 304. Given that support, MOPs can satisfy the requirement of executing progressively at the rate of the inbound dataflow and accomplish efficient Prefetch Pipelining.

As shown, processing core 602 may comprise scanning/indexing PE 622 and XCAM PE 624 as its set of PEs 620. As noted, PEs 620 are the physical entities responsible for executing MOPs, with their underlying UOPs, and for realizing other sophisticated control mechanisms. Various incarnations of processing elements are described herein, where each incarnation supports a distinct subset of the MOP and control space, providing different and distinct functionality from the perspective of query execution. Each of the different PE forms is now addressed where those which support MOPs employing UOP programs implicitly contain a UOP processing engine.

Scanning/Indexing PE 622 implements MOPs which analyze database column groups stored in local memory, performing parallel field extraction and comparison, to generate row pointers (row ids or RIDs) referencing those rows whose value(s) satisfy the applied predicate. For some MOP forms, a metadata Value List (which is an abstract term for a logical tuple list flowing through an IMC) containing a column of potentially sparse row pointers may be given as input, in which case the scan occurs over a sparse subset of the database. For other forms, scanning occurs sequentially over a selected range of rows.

The selection predicate is stipulated through a micro-op (UOP) program of finite length and complexity. For conjunctive predicates which span columns in different column groups, scanning may be done either iteratively or concurrently in dataflow progression through multiple MOPs to produce the final, fully selected row pointer list.

Inasmuch as the Scanning/Indexing PE 622 optimizes scanning parallelism and is capable of constructing and interpreting compacted what are known as bitmap bundles of row pointers (which are a compressed representation of row pointers, sparse or dense, that can be packed into logical tuples flowing through an IMC), it operates most efficiently for highly selective predicates, amplifying the benefits thereof. Regardless, its MOP support locates specific database content.

Scanning/Indexing PE 622 also implements MOPs which project database column groups from HARP memory 304, search and join index structures, and manipulate in-flight metadata flows, composing, merging, reducing, and modifying multi-dimensional lists of intermediate and final results. Depending on the MOP, input is one or more Value Lists whose content may be interpreted in a one- or two-dimensional manner, where two-dimensional lists may have an arbitrary number of columns (which may have arbitrary logical width).

In the context of list reduction, a UOP program of finite length and complexity is stipulated as a predicate function, to qualify one or more components of the input Value List elements, eliminating tuples which do not qualify. List composition involves the combining of related lists into a single output format which explicitly relates the input elements by list locality, while list merging involves intermingling input tuples of like size in an unrelated order. Modification of lists involves a UOP which can generate data-dependent computations, to replace component(s) of each input tuple.

The Scanning/Indexing PE 622 may also be used for joins with indexes, like a Group Index, which involves the association of each input tuple with potentially many related data components, in a one-to-many mapping, as given by referencing the index via a row pointer component contained in each input tuple. MOPs implemented by the Scanning/Indexing PE 622 may thus relate elements of a relational database in by query-specific criteria, which is useful for any query of moderate to advanced complexity.

XCAM PE 624 implements MOPs which perform associative operations, like accumulation and aggregation, sieving, sorting and associative joins. Input is in the form of a two-dimensional metadata Value List which can be interpreted as containing at least two columns related by list locality: key and associated value.

Accumulation occurs over all data of like keys (associatively), applying one of several possible aggregation functions, like Summation or an atomic compare and exchange of the current accumulator value with the input value component. A direct map mode exists which maps the keys directly into HARP memory 304, employing a small cache (not shown) to minimize memory access penalties. A local mode of accumulation exists, as well, to realize zero memory access penalties by opportunistically employing the cache, at the risk of incomplete aggregation.

Sieving involves the progressive capture of keys qualifying as most extreme, according to a programmable sieving function, generating a result list of the original input keys and values such that the last N tuples' keys are the most extreme of all keys in the original input. Iterative application of Sieve can converge on a sorted output, over groups of some small granularity.

Sorting can also be accomplished through construction and traversal of either hashes or B-Trees, which are constructed to relate each input key to its associated value with a structure that is efficient to search and join with.

Within each of PEs 620 thus may be a UOP Processing Engine (not shown). Whereas PEs 620 execute MOPs in a dataflow fashion at the higher levels, embedded UOP Processing Engines in PEs 620 realize the execution of UOPs, which embed within their logical MOP parent to serve its low-level data manipulation and analysis needs. In some embodiments, the UOP processing engine is code-flow logic, where a UOP program is executed repetitively by a parent Processing Element at MOP-imposed boundaries, given MOP-extracted input data, to produce results interpreted by the parent MOP.

Considering the code-flow nature, each UOP engine has its own program storage, persistent register set and execution resources. It is capable, through appropriate UOP instructions, to accept data selected from the parent MOP and to simultaneously execute specified data manipulation or analysis thereon, in combination with some stored register state. In this manner, this tiny code-flow processor is able to fit seamlessly into the dataflow as a variable-latency element which, at the cost of increased latency, is capable of performing any of the most complex low-level data manipulation and analysis functions on the dataflow pouring through. The capability of the MOP to select and present only those data required for UOP processing, at a fine granularity, minimizes the latency imposed by the UOP code flow, maximizing overall dataflow throughput.

FIG. 7—C2 Data Structures

The C2 solution utilizes various hardware-friendly data structures to assist in hardware accelerating database operations by HARP modules 204. As described below in Table 1, the C2 solution may employ three different columns. These columns enable efficient utilization of HARP logic 302 and HARP memory 304, and bandwidth and disk bandwidth with storage infrastructure 112.

TABLE 1 Column Formats Name Description Packed All rows in the column are packed consecutively. All rows Column have a uniform size in each column. A packed column can be disk-resident or memory-resident. The packed Column can be sorted with respect to the primary key. Sorted- This is for the date columns. The date is stored in the Compressed form of running length. Each date has an entry for re- Column cording its starting count and ending count. The offset could be zero if the date does not have any item in the table. It companion column can be accessed with staring and ending counts. Companion The companion table links the Sorted-Compressed Column Column to other columns.

In general, hot columns (i.e., columns having active or frequent access) stay in the HARP memory 304 so that they can be accessed randomly fast. Warm Columns (i.e., columns having less active access) also stay in the HARP memory 304; but occasionally, they may be evicted to a disk in storage infrastructure 112. Cold columns usually be held in storage infrastructure 112, but may be partially brought into HARP memory 304, e.g., for one time usage. In some embodiments, date columns in the Sorted-Compressed format will be held in the memory of host system 202 and accessed by the software running on host 202.

In general, there is a single entry point for HARP module 204 to identify all the database columns. In particular, as shown in FIG. 7, a root table 702 points to all the available table descriptors 704. The table descriptors 704 in turn point to their respective table columns 706. Each table stores multiple columns in the VA memory space. Each of these tables will now be further described.

As noted, root table 702 identifies all the tables accessed by HARP module 204. In some embodiments, each entry in the table takes 8 bytes. When needed, multiple Root Table blocks can be chained by a next pointer. The Descriptor Pointers in the root table 702 points to the individual table descriptors. The indices of the Descriptor Pointers also serve as the table ID. To simplify the hardware design, a CSR (Control Status Register) may be employed to store the Root Table information as long as the hardware accessible Table IDs and Descriptors' information is retained in HARP module 204.

Each database defined table has a table descriptor 704. All the table descriptors 704 may reside in the HARP memory 304. A table descriptor 704 may comprise different groups of data. A group may contain one or more columns. Within a group, the data is organized as rows. A group of data resides in a memory plane which is allocated to it. A data element in a particular plane has direct reference to its corresponding element in another plane. The relationship of the addresses among all the element pairs is the same arithmetical computation. The table descriptor is portable because the present invention utilizes a global virtual address space. In other words, when copying the table descriptor from one virtual memory location to another, all the information in the table is still valid.

In the C2 solution, the data structures of the database are architected to optimize database data processing in HARP hardware. All table columns/column groups, indices and meta-data are defined in a global database virtual address space (DBVA). A reserved DBVA section is allocated for table descriptors 704 as part of the meta-data. Table descriptors 704 include information about a table, such as the table name, number of rows, number of columns/column groups, column names, width(s) within a column group, etc. In addition to the information of data layout and access information in the VA space, the table descriptors 704 also have information about the compression types/algorithms used for each individual column. In the present invention, hardware can directly use this information to accomplish database queries and table element insertion, update, and deletion.

FIG. 8—Table Column Layout

FIG. 8 is now provided to provide further detail on the structure of a table in column-store format as employed by the C2 solution of the present invention. As shown, each database table is broken into multiple columns or column groups having a fixed width. Variable width columns are also supported using column hopping or a column heap structure with linked lists. In the C2 solution, a column group can have one or more columns packed together. Because of the simple arithmetic mapping or the single indirection in the companion column, the hardware and software of the present invention can easily access rows across the columns without any degradation in performance; thus, the C2 solution can provide the same functionality and benefits as known row store databases. Table and column descriptors may also be embedded in the MOPs and query tasks.

Of note, in the present invention, the columns or column groups possess an implicit row id (RID). A RID is considered implicit because it is not materialized as a part of a column or column group. Instead, each column and column group is designated a starting RID, which corresponds to an address in the global database virtual address space, which is then mapped to a physical address in HARP memory 304. Since each column and column group is a fixed width, the RID can provide the basis for arithmetically calculating the memory address of any data in the column or column group.

In some embodiments, all columns are packed together in the single DBVA. In addition, a meta-data structure may be employed to facilitate certain column accesses. For example, as shown, a row pointer primary key index may comprise a sorted list of primary keys and their associated row id (RID) in a column or column group. Of course, a B-tree index may be used as an alternative to this type of index.

In the present invention, two active sets of database regions are maintained, i.e., a main database region and an augment region for newly added data. Query processing operates on both regions and is accelerated by the HARP module 204. The augment region is utilized to hold new inserted items. Optionally, the augment region may be rolled into the main region. For example, as shown in FIG. 8, RIDs 1−n are the main region, while RIDs n+1, etc. comprise the augment region.

Deletion updates may be committed into the main region right away. To alleviate the drastic changes across all the columns in a table, the present invention may allocate a valid or invalid bit. A row deletion in a table, therefore, becomes a trivial task of setting the appropriate bit in every column group in the table.

FIG. 9—Example of a SQL Query

FIG. 9 shows one of the 22 TPC-H queries, query #3, and how it would be executed using the machine code database instructions. TPC-H queries are published by the Transaction Processing Performance Council (TPC), which is a non-profit organization to define benchmarks and to disseminate objective, verifiable TPC performance data to the industry. TPC benchmarks are widely used today in evaluating the performance of computer systems. This particular query is a shipping priority query to find the potential revenue and shipping priority of the orders having the largest revenue among those that had not been shipped of a given date. The market segment and date are randomly generated from the prescribed range, and “BUILDING” and Mar. 15, 1995 are the example here. This query is a complex multiple table join of three tables, CUSTOMER, ORDERS, and LINEITEM tables.

C2 Software 110 will decompose this query into 24 MOPs to send to HARP module 204, along with their dependency information, which establishes the topology of the dataflow from MOP to MOP. All MOPs are started and hardware processing begins in pipelined fashion, with each MOP's results being fed to one or more downstream consumers over one or more dedicated logical IMC connections.

The responsibility of the first MOP, ScanCol(0), is to reference HARP memory 304 to find all the customers in the CUSTOMER table who belong to the ‘BUILDING’ market segment, producing into IMCO all matching CUSTOMER references in the form of one RID per qualified row. Revindex(1) then traverses a reverse index residing in 304, pre-built to relate customers to their one or more orders residing in the ORDERS table, outputting references to all orders made by the given customers. Because the CUSTOMER references are no longer necessary and to boost performance by reducing utilization of IMC transmission resources over IMC2, the ListProject(2) removes the original customer references after the reverse index join, leaving only the ORDER references. The ScanRPL(3) MOP then scans these orders' 0 ORDERDATE column, retaining ORDER references only to those orders whose order date occurs before the date ‘1995-03-15’.

Progressing onward through IMC3, the dataflow entering Revindex(4) consists of ORDER table references (RIDs) which have satisfied all criteria mentioned thus far: each order was placed by a customer in the ‘BUILDING’ market segment before the date Mar. 15, 1995. To finish evaluating the “WHERE” clause of the illustrated SQL query statement, these orders must be qualified in terms of certain properties of their related line items.

The purpose of the Revindex(4) MOP is then to associate each of the qualifying orders to its one or more constituent line items from the LINEITEM table, returning appropriate references thereto. At this point, the flow contains a two-column tuple list relating ORDER references (RIDs) to LINEITEM RIDs, multicasting identical copies of these tuples into IMC4 and IMC5. ListProject(5) extracts only the LINEITEM RID column from the dataflow in preparation for ProjRpl(6), which extracts each line item's L_SHIPDATE column value, feeding these ship dates to IMC7. ListCompose(7) consumes IMC7 along with IMC5, executing a composition of the input lists to create a three-column tuple list where each tuple contains an ORDER RID, an associated LINEITEM RID and its ship date. ListSelect(8) consumes the composed list from IMC 8 and selects only those tuples having ship date older than ‘1995-03-15’, thus completing the “WHERE” clause requirements.

Again, at the output of ListSelect(8), the dataflow still logically appears as a three-column tuple list where each tuple relates an ORDER RID to one of its associated LINEITEM RIDs and that line item's ship date. It should be noted in this flow that multiple distinct LINEITEM RIDs may appear (in different tuples) with an identical ORDER RID, a definite possibility here since a single order may be comprised of an arbitrary number of line items in the target database and this query specifically requests only those line items satisfying the ship date criteria. The redundancy of ORDER RIDs in the list suggests an aggregation step will be needed to realize the SUM of the SQL select statement, but before that, some more data must be gathered and calculations done.

IMC9 and IMC10 both carry the output of ListSelect(8), identically. ListProject(9) extracts only the LINEITEM RID column from IMC9, passing that on to both ProjRpl(12) and ProjRpl(11), which fetch each referenced LINEITEM's L_EXTENDEDPRICE and L DISCOUNT, respectively. Those procured extended price and discount data are then composed together by ListCompose(13) to form a two-column tuple to be carried via IMC17. ListTupleArith(14) implements the arithmetic process of computing (L_EXTENDEDPRICE*(1−L_DISCOUNT)) on a per-tuple basis before sending this arithmetic result to ListCompose(15). In the meantime, ListProject(10) extracts the ORDER RID column from the output of ListSelect(8), such that ListCompose(15) can make a two-column composition relating, within each tuple, an ORDER RID to its line item's arithmetic product.

The final hardware step to complete the query involves fully evaluating the SELECT clause, including its SUM aggregation function. The remainder of the MOP flow of FIG. 9, beginning with the output of ListCompose(15), is dedicated to this process.

AssocAccumSum(16) receives from IMC19 with each of the two-column tuples relating an ORDER RID to one of its line item's (L_EXTENDEDPRICE*(1−L DISCOUNT))product, computing a summation of these values independently for each distinct ORDER RID. For example, a given ORDER RID may appear twice in IMC19 (once in two different tuples), having two distinct LINEITEMs which satisfied all criteria thus far. Each of these LINEITEMs would have generated its own product in ListTupleArith(14), such that the aggregation process of AssocAccumSum(16) must sum them together. The result is a distinct sum of products over each distinct ORDER RID, realizing the SQL SUM aggregation function, here named REVENUE within the query.

Once the aggregation has completed for a given ORDER RID, ListProject(17) extracts the ORDER RID itself, passing it to ProjRpl(18), ProjRpl(19) and ProjRpl(20). These MOPs gather in parallel the referenced orders' 0 ORDERDATE, O_SHIPPRIORITY, and 0 ORDERKEY, repectively, while ListCompose(21) forms a two-column tuple consisting of 0_SHIPPRIORITY and 0 ORDERKEY. ListCompose(22) meanwhile forms a two-column tuple comprised of O_ORDERKEY and REVENUE. The final MOP, ListCompose(23), composes the two two-column tuple lists into a final four-column tuple list which satisfies the SQL query and its SELECT statement.

It should be noted in this example that the SQL query SELECT actually stipulates L_ORDERKEY. But an optimization may be applied here, knowing that O_ORDERKEY is functionally equivalent when used in this manner, thus avoiding the need to carry any LINEITEM RIDs beyond IMC11 or IMC12.

FIG. 9—Example of a Dataflow through the HARP

In FIG. 9 we have described how an SQL statement gets mapped into a logical MOP DAG (directed acyclic graph) which gets executed in a dataflow fashion with IMC chaining between MOPs. FIG. 10 illustrates an exemplary dataflow through PEs 620 in HARP logic 302 for the same TPC-H SQL #3 query shown in FIG. 9. As noted, C2 Software 110 will decompose this query task into 10 PE stages to send to HARP module 204, along with their MOP and UOP instructions and dependency information.

Stage 1 is performed by Scanning PE 1002 is to find all the customers in CUSTOMER table that is in “BUILDING” market segment and passes the results (C_RIDs of matching customer records) in an IMC to Indexing PE 1004.

Stage 2 is a join operation of C_CUSTKEY=O_CUSTKEY performed by Indexing PE 1004 using a reverse index method. Each C_RID of Stage 1's matching customer records corresponds to an O_RID hitlist of ORDER table records, given a customer may place multiple orders. The results (O_RIDs) are passed in an IMC to Scanning PE 1006.

Stage 3 is performed by Scanning PE 1006 to read the 0 ORDERDATE field of all the matching orders (ORIDs) that Stage 2 outputs, compare for “<‘1995-03-15”, and passes the results (O_RIDs) in an IMC to Indexing PE 1008.

Stage 4 is a join operation of O_ORDERKEY=LORDERKEY performed by Indexing PE 1008 using a reverse index method. Each O_RID of Stage 3's matching order records corresponds to an L_RID hitlist of LINEITEM table records, given an order may have multiple line items. The results (L_RIDs) are passed in an IMC to Scanning PE 1010.

Stage 5 is performed by Scanning PE 1010 to read the L_SHIPDATE field of all matching line items (L_RIDs) that Stage 4 outputs, compare for “>‘1995-03-15”, and passes the results (L_RIDs) in 3 IMCs to Indexing PE 1012, 1014, and 1016.

Stage 6 is a column extraction/projection operation done by Indexing PE 1012, 1014, and 1016 to get L_ORDERKEY, L_EXTENDEDPRICE, and L_DISCOUNT column.

Stage 7 is a list merge operation of 2 columns (L_EXTENDEDPRICE and L_DISCOUNT) done by Indexing PE 1018.

Stage 8 is an aggregation operation of REVENUE of each L_ORDERKEY group, done by XCAM PE 1020 based on outputs of Indexing PE 1012 and 1018. As the SQL statement defines, REVENUE is calculated as the sum of (L_EXTENDEDPRICE*(1-L_DISCOUNT)). Note that even though the GROUP BY defines the group key as concatenation of L_ORDERKEY, O_ORDERDATE, O_SHIPPRIORITY, the group key is simplified to L_ORDERKEY since it is already a unique identifier. The output of XCAM PE 1020 is a pair list of group key (L_ORDERKEY) with its REVENUE.

Stage 9, done by Indexing PE 1022 and 1024, is a column extraction of O_ORDERDATE based on L_ORDERKEY output of XCAM PE 1020.

Stage 10, done by XCAM PE 1026, is a sieve (ORDER BY) operation of REVENUE, O_ORDERDATE to output top N groups with largest REVENUEs. These outputs are placed at a “result” buffer area in HARP memory 304, ready to be retrieved by DBMS software 114.

FIG. 10—Example of a Dataflow through the HARP

FIG. 10 illustrates an exemplary dataflow through PEs 620 in HARP logic 302 for the same TPC-H SQL #3 query shown in FIG. 9. As noted, C2 Software 110 will decompose this query into 7 query tasks to send to HARP module 204, along with their task command and dependency information.

Task 1 is performed by a Scanning PE is to find all the customers in CUSTOMER table that is in BUILDING market segment and passes the results in an IMC to an Indexing PE.

Task 2 is performed by an Indexing PE to extract the C_CUSTKEY field (column) of Task 1's matching customer records and search against the 0 CUSTKEY index table and passes the results in an IMC to a Scanning PE.

Task 3 is performed by the Scanning PE to read the O_ORDERDATE field of all the matching orders that Task 2 returns and passes the results in an IMC. Given a customer could have placed multiple orders, Task 2's CUSTKEY search could result in multi-record hit lists. An Indexing PE then extracts the appropriate records and passes the results in an IMC to a Scanning PE.

Task 4 is performed by a Scanning PE to compare the 0 ORDERDATE of each order record from Task 3 and returns only those records with dates before Mar. 15, 1995 in an IMC to a set of Indexing PEs.

Task 5 is performed by a set of Indexing PEs and then an XCAM PE to extract the 0 ORDERKEY field of Task 4's qualified orders and search against the L_ORDERKEY index table.

Task 6 is performed by a set of Indexing PEs to extract the L SHIPDATE field of all the matching line items that Task 5 returns.

Task 7 is then performed by an XCAM PE to compare the L SHIPDATE of each line item record from Task 6 and sort those records with shipping date later than Mar. 15, 1995. The sorted line item records are placed at a result buffer area in HARP memory 304, ready to be retrieved by DBMS software 114.

Explanation of the Global Database Virtual Addressing Scheme.

As noted, the C2 solution of the present invention can employ an arbitrarily large virtual address space (DBVA) to globally address an entire database. This addressing scheme is utilized to assist hardware acceleration, because it allows HARP module 204 and C2 software 110 to arithmetically calculate the location of any data in a database. FIG. 11 illustrates an exemplary logic flow of the addressing schemes employed and FIG. 12 illustrates a structure of the global database virtual address scheme of the present invention.

Referring now to FIG. 11, when a query is received from MySQL software 114, it is eventually submitted to query execution engine 412 for processing. At this level, execution engine 412 operates on the database using what are known as logical addresses. A logical address is utilizes the naming convention of the SQL query. For example, a logical address may consist of a database name, a table name, a column group name, and a column name which is then converted to a DBVA. Since it is a virtual address, the DBVA may be arbitrarily large, e.g., in excess of a petabyte.

The DBVA may then be converted to physical addresses of the data in storage 112 or in the memory of host system 202. For example, as shown, a storage manager 420 may have a mapping that converts a DBVA to a storage physical address (SPA). This allows for retrieval of data from storage 112. In addition, buffer manager 414 may convert the DBVA to a physical address in the memory of host system 202, for example, using a standard memory manager in the operating system running on host system 202. Such conversions are well known to those skilled in the art.

However, of note, the present invention may employ an extensible addressing scheme for one or more HARP modules 204. As noted, HARP modules 204 may comprise a relatively large memory, i.e., in excess of 256 gigabytes or more. In addition, the C2 solution may comprise nodes that have multiple HARP modules 204 and may also support nodes. Accordingly, the present invention may comprise a novel, layered virtual addressing scheme.

In one embodiment, each HARP module 204 translates the DBVA to a physical address (HPA) of HARP memory 304. For example, HARP module 204 may utilize a translation-lookaside buffer (TLB) for this translation.

Alternatively, especially in a multi-HARP module environment or federated system, each HARP module 204 may employ a secondary virtual address space (called an HVA) that is underlies the DBVA. Referring now to FIG. 12, in this embodiment, memory manager 418 may thus include a DBVA-HVA service that maps a DBVA address into yet another virtual address in the HVA. Memory manager 418 may then translate the HVA into a HPA using, for example, a TLB for the respective HARP memory 304.

This feature allows the present invention to continue utilizing a single global virtual address space for a database even where system 100 includes multiple HARP modules 204 or multiple nodes of HARP modules. Accordingly, systems of the C2 solution can continue to employ hardware-friendly addressing for hardware acceleration of the database operations regardless of the number of HARP modules 204 or the size of the database.

Multi-Versioning and Continuous Updates

Embodiments of the present invention can provide fine grain concurrency control in the presence of database updates. For read-only transactions, historical snapshots of the data in pages in HARP memory 304 are created. In addition, historical snapshots can be created for pages in the memory of host 202 and in storage 112. For illustrative purposes, the following description is directed towards multi-versioning and continuous updates of data in various pages of HARP memory 304. However, one skilled in the art will recognize how the concepts of the present invention for multi-versioning and continuous updates can be extended to host 202 and storage 112.

The following description begins by generally explaining various concepts of multi-versioning and continuous updates that may be implemented by embodiments of the present invention. FIGS. 11 and 13 illustrate an example of how multi-versioning and continuous updates may be accomplished.

Embodiments of the present invention provide fine grain concurrency control for transactions in the presence of database updates. During operations, each transaction is assigned a snapshot version number or SVN. A SVN refers to a historical snapshot of the database that can be created periodically or on demand. Transactions are thus tied to a particular SVN, such as, when the transaction was created. Queries belonging to the transactions can access data that is consistent as of a point in time, for example, corresponding to the latest SVN when the transaction was created.

At various times, data from the database stored in a memory can be updated using the snapshot data corresponding to a SVN. For example, when a transaction is committed, a snapshot of the database with a new SVN is created based on the data modified by the transaction and the snapshot is synchronized to the memory. As another example, when a transaction query requires data from a version of the database corresponding to a SVN, the data in the memory may be synchronized with the snapshot data corresponding to that SVN. This feature, in essence, creates multiple versions of the database and allows read-intensive database applications to operate in the presence of database updates.

A query can be isolated to a consistent historical snapshot. Changes made by committed update transactions may be visible to other transactions unless they are operating on a historical snapshot. Queries and data manipulation are executed within a transaction boundary. In other words, a transaction serves to provide a consistent read for its related queries and query tasks. When multiple queries are executed within the same transaction, these queries will also operate on the same version. Two modes of transactions, serializable and snapshot serializable may be supported.

A serializable mode transaction is exposed to the latest committed changes. Snapshot serializable transactions are exposed only to the snapshot version they were assigned when they started. If a transaction does not exist when the query/data manipulation is executed, C2 software 110 may start a transaction automatically and auto commit when the statement is processed.

For locking, C2 software 110 may rely on MySQL software 114 for table locking. Internally, memory manager 418 can support page level locking in HARP memory 304 which can be used to provide more concurrency if required. For serializable transactions, write locks are issued so that no other serializable transaction can access the same data. For snapshot serializable transactions, transactions are not blocked on account of the serializable transactions.

In the present invention, there are several embodiments for handling updates during the data manipulation (DML) processing. In one approach, changes made during the DML processing are not made to the data inline, but stored in offline auxiliary structures. This allows queries to run interrupted.

Alternatively, a second approach may be employed that involves keeping shadow copies of pages touched by transactions, but maintaining a different version until the changes are synced. If the pages need to be flushed to storage 112 after changing in the memory of host system 202, the storage manager 420 may maintain two versions of the pages, one for the system SVN and the next version. In this approach, the queries are relatively uninterrupted since the version needed by them is not affected by the transactions.

Another approach can utilize a change log. In some embodiments, the change log can be in the form of undo or redo operations. In this approach, the changes are made directly to the pages in the buffer cache in the same page (rather than a shadow copy). These changes are made in the buffer cache and may get flushed to storage 112 but are not made in HARP memory 304 until prefetched for query processing. A status map may be kept by C2 software 110 to indicate the status of data on memory in host system 202, storage 112, and HARP memory 304. Thus, when a prefetch request is made for certain data, this map is consulted to ensure that data is present in the version required and if not, a change log can be used to construct it.

Non-Blocking Continuous Updates by the C2 Solution

Referring now to FIGS. 13-14, an embodiment of continuous updating by the C2 solution is conceptually illustrated. As shown in FIG. 13, a timeline of query transactions X, Y, and Z are shown. As time progresses, snapshots are generated periodically (or at commit time for sync on commit).

Snapshots may be managed by a snapshot manager in utilities 426 and, as noted, are identified by their SVN. In some embodiments, the snapshots may be managed by update manager 424. Snapshots can be created in various ways. In snapshot on commit, a snapshot is created from the data modified by a transaction at commit time. In a timed snapshot, several transactions' committed data is buffered and associated with a snapshot based on a timer, which can be user specified.

A serializable mode transaction does not bind to an SVN. A snapshot mode transaction binds to an SVN. Pages are associated with a specific SVN. A newly created snapshot transaction is associated with the current SVN. A query task inherits the SVN from the transaction it belongs to. All the tasks belonging to a query will inherit that SVN too.

Synching is the process of bringing a page to the required SVN. Pages can be on different SVNs. In order for a task to execute, it's required that all the pages touched by that task are present at the correct SVN. For example, as shown, for various pages in HARP memory 304, update manager 424 manages a snapshot version number (SVN), e.g., SVN2, SVN3, and SVN4 and a set of undo and redo tables for each SVN. Each transaction is associated with the SVN of the current snapshot at the time it is created. This SVN is used for all queries and tasks executed for that transaction. Snapshots provide approximate up-to-date data for queries while preserving performance since many queries and tasks may work on the same snapshot of data. In some embodiments, the snapshot data structures may be stored in B+ trees. In these data structures, uncommitted data of active transactions and committed data since the last snapshot are tracked.

In some embodiments, the SVN may simply be a value that is monotonically increasing and assigned by update manager 424. SVNs can be recycled when the system restarts. In some embodiments, there is a single global snapshot version number maintained by update manager 424 that all snapshot serializable queries can access. Serializable transactions see the version of data for a page in HARP memory 204 that includes committed changes that have occurred since the last sync. When the committed changes are synced, the SVN is incremented.

When a transaction starts, it is assigned a snapshot version (the current system SVN). This is the version of the data that all queries within this transaction will see for the duration of the transaction. Hence, as shown in FIG. 13, query transaction X is assigned SVN1 since it started during SVN1. Likewise, query transactions Y and Z are assigned SVN3 and SVN4, respectively. Of note, as shown, query transactions may span across multiple SVNs, but may be referenced according to the SVN in which they started.

A query executes on the snapshot with the transaction's SVN. For example, as shown in FIG. 13, the three shown query tasks in query transaction X operate on data from SVN 1. The query tasks in transaction Y operate on data from SVN2, and query tasks in transaction Z will operate on data from SVN4. The transaction's SVN is determined at transaction start time. If a transaction does not exist then one may be automatically created that lasts for the duration of the query.

However, as shown, some tasks may be submitted at a different SVN from the SVN assigned to their transaction. For example, transaction X may include a task that starts in SVN2 rather than SVN1. Transaction Y may include a task that starts in SVN3 rather than SVN2. These tasks may thus need to have one or more pages synched to a different SVN, which will now be explained with reference to FIGS. 14-16.

Referring now to FIG. 14, transitions between SVNs may be mapped using redo and undo tables that are logged in the snapshot data structures. As shown, a page from HARP memory 304 may be synched between SVN2 and SVN3, SVN3 and SVN4, and SVN2 and SVN4. The redo table is used to convert the data in an older snapshot to a newer snapshot. The undo table is used to convert the data from a newer snapshot to an older snapshot. The redo table from a snapshot can also be maintained to the next snapshot and the undo table from a given snapshot to the prior snapshot. This data thus allows for reconstruction of snapshots with minimal overhead.

In some embodiments, these tables are implemented in 32 byte chunks. These tables may map a VA to a 32 byte chunk of data. The VA is a 32 byte-aligned address. The VAs may be in sorted order.

The task manager will execute a prefetch task to load the data required by the query or task for its SVN. The prefetch task is submitted as part of the query execution plan generated by query plan generator 410. The memory manager 418 maintains page to SVN mapping. If a page that needs to be prefetched is missing or at the wrong SVN, then the page may be updated to the correct SVN.

During operation, tasks for queries that are part of a transaction are submitted for execution on HARP module 204. These tasks are submitted with the SVN required for that task so that the task manager 416 can handle the task appropriately when it reorders this task across sync barriers. In addition, elsewhere, queries and query tasks in these transactions may request blocks of this SVN from the memory in host system 202.

A snapshot serializable transaction starting at SVN2 will be shown those HARP memory pages having a SVN2. Meanwhile, other transactions may make changes to data after this point and the pages touched by these transactions will be marked with a different SVN, such as SVN3 or SVN4, when they commit.

A transaction starting in serializable mode may be able to see these changes that have not been synced yet. Since all committed changes for multiple transactions go into the same SVN, a transaction in serializable mode when it starts must see a consistent read throughout its lifetime. For example, if a query task starts in serializable mode and can see committed changes which are part of SVN2, another query task may come along and make more changes into SVN2 but these will not be visible to that first query task.

In the present invention, a way to solve this is for query tasks to acquire its table level lock in write mode even if it just wants to do queries. This does not impact snapshot serializable queries which only lock in read mode and can always progress. Alternately, the locks held by memory manager 418 can be held for transaction duration which will achieve the same effect.

Then pages are being changed, their state across the different components may need to be tracked. As noted, in some embodiments, C2 software maintains track of the SVNs of pages in storage 112, host system 202, and HARP memory 204. In some embodiments, a page map is implemented in the form of a hash table having the address prefix of each page.

In some embodiments, an isolation level may specify that all transactions occur in a completely isolated fashion; i.e., as if all transactions in the system had executed serially, one after the other. Read committed isolation can be supported by releasing a read lock on a table at the completion of the read.

FIG. 15 illustrates an example of converting a page of HARP memory 304 from SVN2 to SVN3 based on information in a redo table. As shown, HARP logic 302 may utilize one of its XCAM PEs 624 to transition a page of HARP memory 304 between SVN2 to SVN3. In particular, a redo table may be fed into XCAM PE 624, which then performs an associative accumulate function to change values in the page of HARP memory 304. Further details of this function are shown with reference to FIGS. 17 and 18.

FIG. 16 illustrates an example of converting a page of HARP memory 304 from SVN3 to SVN2 based on information in an undo table. As shown, HARP logic 302 may utilize one of its XCAM PEs 624 to transition a page of HARP memory 304 between SVN3 to SVN2. In particular, an undo table may be fed into XCAM PE 624, which then performs an associative accumulate function to change values in the page of HARP memory 304. Further details of this function are shown with reference to FIGS. 17 and 18.

FIGS. 17-18 illustrate associative accumulation machine code database instructions that may be employed to convert pages between SVNs. In particular, FIG. 17 illustrates a full replace function, while FIG. 18 illustrates a selective replace function.

In general, an associative operation implemented by XCAM PE 624 involves the writing of each input value, mapped by its associated key value, to a corresponding location in HARP memory 304. Since the input keys may vary arbitrarily in their order and value, this process can implement a scattering of writes to desired portions of HARP memory 304. The scattering writes may be performed in a manner particularly useful for making anything from large, sweeping to small, incremental updates of database column contents. As noted, at least two approaches may be available: full replace shown in FIG. 17 and selective replace shown in FIG. 18.

The full replace function (or coalescing form) shown in FIG. 17 first collects writes to adjacent memory locations over a window of address space, before dispatching them to HARP memory 304. While avoiding reads of original memory contents, coalescing takes advantage of the potential for spatial locality in the input key stream, thereby minimizing the write and nullifying the read bandwidth required, to provide a high performance solution for executing updates of memory 304. Since reads are avoided and since the granularity of the update size can be much less than that of the coalescing window, any key location not addressed within the window may have its associated memory contents cleared to a constant value or assigned otherwise undefined data. A small cache may be utilized to perform the coalescing.

Referring now to FIG. 18, the selective replace function is shown. This function may also be considered a non-coalescing form and may be more precise and selective in updating HARP memory 304, preserving data in adjacent locations, but may utilize higher memory read and write bandwidth requirements. This approach is especially useful for small update granularity (e.g. 1-bit) where spatial locality cannot be guaranteed among the updates requested in the input stream and preservation of surrounding data is necessary. Again, a small cache may be used to mitigate unnecessary memory accesses, in an opportunistic manner.

Other embodiments of the invention will be apparent to those skilled in the art from consideration of the specification and practice of the invention disclosed herein. It is intended that the specification and examples be considered as exemplary only, with a true scope and spirit of the invention being indicated by the following claims. 

What is claimed is:
 1. A method of versioning data in a database maintained within a host computer system coupled with a hardware accelerator, the host computer system including a host memory, and the hardware accelerator including a hardware accelerator memory, said method comprising: determining snapshots of data in the database at an interval; identifying differences in the data; determining a set of changes that takes the data from one snapshot to another snapshot; and generating a mask comprising the set of changes that is applied by a machine code database instruction for modifying the data in pages of the hardware accelerator memory from one snapshot to another snapshot.
 2. The method of claim 1, wherein generating the mask comprises generating a mask having a fixed width format that is independent of the data type being modified.
 3. The method of claim 1, wherein generating the mask comprises generating a mask having data that is in a same compressed form as the data in the database.
 4. The method of claim 1, further comprising storing a mask for an active snapshot in the host memory and storing at least one additional mask for completed snapshots in the host memory.
 5. The method of claim 1, further comprising storing a mask for an active snapshot in the host memory and storing at least one additional mask for completed snapshots in the hardware accelerator memory.
 6. The method of claim 1, further comprising setting the interval based on a configuration.
 7. The method of claim 1, further comprising setting the interval based on a size of accumulated changes.
 8. The method of claim 1, further comprising applying the mask to data in the hardware accelerator memory upon a transaction commit.
 9. The method of claim 1, further comprising applying the mask to data in the hardware accelerator memory upon demand by a query.
 10. A tangible, non-transitory, computer readable medium comprising program code for performing a method of versioning data in a database maintained within a host computer system coupled with a hardware accelerator, the host computer system including a host memory, and the hardware accelerator including a hardware accelerator memory, said medium comprising: program code for determining snapshots of data in the database at an interval; program code for identifying differences in the data; program code for determining a set of changes that takes the data from one snapshot to another snapshot; and program code for generating a mask comprising the set of changes that is applied by a machine code database instruction for modifying the data in pages of the hardware accelerator memory from one snapshot to another snapshot.
 11. The computer readable medium of claim 10, wherein the program code for generating the mask comprises program code for generating a mask having a fixed width format that is independent of the data type being modified.
 12. The computer readable medium of claim 10, wherein the program code for generating the mask comprises program code for generating a mask having data that is in a same compressed form as the data in the database.
 13. The computer readable medium of claim 10, further comprising program code for storing a mask for an active snapshot in the host memory and program code for storing at least one additional mask for completed snapshots in the host memory.
 14. The computer readable medium of claim 10, further comprising program code for storing a mask for an active snapshot in the host memory and program code for storing at least one additional mask for completed snapshots in the hardware accelerator memory.
 15. The computer readable medium of claim 10, further comprising program code for setting the interval based on a configuration.
 16. The computer readable medium of claim 10, further comprising program code for setting the interval based on a size of accumulated changes.
 17. The computer readable medium of claim 10, further comprising program code for applying the mask to data in the hardware accelerator memory upon a transaction commit.
 18. The computer readable medium of claim 10, further comprising program code for applying the mask to data in the hardware accelerator upon demand by a query. 